Systems and methods for a cache-sensitive index using partial keys

ABSTRACT

Systems and methods are disclosed for a cache-sensitive index that uses fixed-size partial keys. The index may include a node comprising a child group pointer, a number of partial keys and a similar number of full-key pointers. The node may also include a record count. The nodes are organized into groups. The groups may contain a number of nodes one greater than the number of partial keys in a node and the nodes in a group may be stored contiguously in memory. The child group pointer and the number of partial keys may fit within a cache line. A method is disclosed for traversing the index, for bulk-loading the index, and for live deletion of records from the index.

FIELD

This disclosure is generally directed to systems and methods forcreating an index supporting fast main-memory database lookups and, moreparticularly, to systems and methods for creating a cache-sensitiveindex that also uses partial keys.

BACKGROUND

Relational databases store data using tables. Generally, a table in arelational database consists of data organized using columns and rows.The columns represent a particular field, such as “last name” or “ID.”The rows represent data records stored in the columns, such as “Smith”and “12345.” A particular table may have millions of rows of data,making a search for a particular row slow and cumbersome. To speedaccess to the records in a table, databases index the rows with an indexstructure having algorithmic search properties. Historically, relationaldatabases have used an index structure, called a B+ tree, to provide theshortest path possible to the desired data. In a B+ tree, a search isperformed from the root of the tree through intermediate nodes down toleaf nodes. The root node and the intermediate nodes are collectivelyknown as index nodes and point to other index nodes or to the leafnodes. The leaf nodes point directly to records in the database (the rowdata). A B+ tree contains copies of the keys in the nodes and has a highnumber of children per node, making the path from the root node to theleaf nodes short. A short path is desirable because it results in feweraccesses to a disk storing the index. Disk accesses have a much sloweraccess time than main memory accesses, but because of the cost of mainmemory, the tables and indexes are generally stored on disk-type storagedevices.

As main memory decreased in price databases stored in main memory becamepractical. Because disk accesses are not a concern of main memorydatabases, the index for a main memory database would optimally seek tooptimize cache memory usage rather than reduce disk accesses. Toaccommodate this, two techniques arose to increase the searchperformance of the B+ tree. The first decreases the number of “down”pointers in a node by addressing groups of nodes rather than individualchild nodes. In such an index, the child nodes are stored contiguouslyin main memory in a node group. Thus, only one pointer is needed topoint to the child nodes even though the node group contains severalchildren. The location of each child (index key) can be calculated bysimple arithmetic, since the nodes contain a fixed number of bytes andare contiguous. The node groups are also organized to avoid crossingcache-lines (e.g. if a particular cache is organized in 64 byte blocks,the cache line falls between every 64 bytes). This technique allowsnodes to contain more key pointers, increasing processor cacheefficiency. The root group of such a cache-sensitive index contains asingle node, but subsequent groups may have one node more than the keysin the parent group. A cache-sensitive index focuses on reducing pointeroverhead (i.e. reducing the number of pointers) and improving spaceutilization so that more keys can be added to the same-sized node. Thistrades off search speed for update speed because updates involve copyingentire groups of nodes rather than individual nodes. FIG. 1 depicts acache-sensitive index with node groups represented by a dashedrectangle.

The second cache-conscious version of the B+ tree is a partial keyindex. A partial key index reduces the size of the index by only storingpartial keys and not full keys in the index nodes. Each node contains aset of down pointers, which point to other nodes or to records (rows).The nodes also contain partial keys and pointers to the full key, whichis located in the record itself. The partial key information includes atwo-byte offset indicating at what position the partial key differs fromthe base key, and two bytes of data after the offset that differ fromthe previous key. For example, if a base key contains “ABCDEF” and thenext key contains “ABEGXY”, the partial key contains an offset of 2 andthe 2 bytes of differing data contain “EG.” Thus, the partial key onlycontains the position of the key that differs from the base and the twobytes of data that differ from the base. A partial key index focuses onlowering key-comparison cost rather than reducing pointer overhead. FIG.2 represents a partial-key index.

The Domain Name System (DNS) uses a distributed network of name servers(lookup nodes) to translate text-based web addresses, such as“www.acme-co.com,” to Internet protocol (IP) addresses, such as“234.562.55.3.” When an Internet user requests a web address, one ormore name servers process the DNS request by looking up the web addressin a database of registered domains. When the name server locates theweb address in the database, the IP address is sent back to the user'scomputing device.

Some name servers must handle millions of DNS requests each second.Furthermore, the name servers must perform the resolution quickly toenhance the user experience on the Internet. Therefore, name servers mayuse main-memory databases to store the records needed to successfullyresolve a DNS request to allow faster access to the data. Furthermore,web addresses are added to and removed from the name server databasedaily. To accurately resolve a DNS request, the name server must rely onan index updated in real time.

Because the traditional DNS resolution process is vulnerable to hacking(i.e. forged DNS data), the industry has begun to implement a secureversion of DNS named DNSSEC (DNS Security Extensions). DNSSEC requireseach DNS lookup node to authenticate the DNS request, thus ensuring thatthe request will not be misdirected to a fraudulent site.

To authenticate a DNSSEC request, the lookup node must determine wherethe web address falls in relation to the DNS zone. Internet addressesare divided into DNS zones in a hierarchical tree-like fashion. The rootzone includes all top-level international, ISO country-code, and genericdomains and are serviced by root name servers. Below the root zone aretop-level domains (TLDs), such as “.com,” “.net,” and “.org.” The TLDsmay be further divided into zones managed by organizations that registerthe second-level domains. These organizations may decide to delegateauthority for sub-zones within lower-level domains. Thus, there may beseveral name servers responsible for the different zones associated witha web address. DNSSEC requires a name server to determine not only thata particular web address exists, but also what falls just after it andprior to it in the zone.

Therefore, it is desirable to introduce an index structure thatfacilitates faster access to large main-memory databases while stillretaining the ability to add and delete records from the index in realtime.

SUMMARY

Disclosed embodiments provide a cache-sensitive index that also containspartial keys, as well as methods for maintaining such an index inreal-time so that the index can be used for lookups that require theindex to reflect live updates to the database. The index combines acache-sensitive index and a partial key index to improve lookup speed.Each node of the index contains a pointer to a sub-group, fixed-sizepartial keys, and full key pointers. The node is configured so that thefixed-size partial keys fit in the first cache line. The full keypointers and key count take up the remaining space in the node, suchthat a node fits in two processor cache-lines.

An ordered index, such as the cache-sensitive index using partial keysdescribed herein, is especially valuable for processes that require notonly a lookup of a record in a database, but also a lookup of the “next”and possibly the “previous” entries once the record is found. Forexample, a DNSSEC proof may require that the DNS server locate not onlya requested domain but also the “next” and “previous” domains sortedaccording to domain naming rules. An ordered index, such as thecache-sensitive index using partial keys, simplifies and expedites thistype of operation by storing ordered index records contiguously inmemory. Because a DNS server processes millions of requests per second,the cache-sensitive index using partial-keys dramatically increasesresponse time for DNSSEC requests.

Consistent with disclosed embodiments, an index for a database isprovided that comprises groups of nodes. The nodes of each group may bestored contiguously in memory. The nodes may comprise a child grouppointer, a number of fixed size partial keys, and a number of full keypointers corresponding to the number of fixed size partial keys. Thenumber of nodes in each group may be one higher than the number of fixedsize partial keys in each node. The partial keys may fit in a cacheline.

Consistent with disclosed embodiments, a method is provided for locatinga search value in a database using such an index. The method may includecomputing a partial key for the search value and starting at the rootnode by setting a current node to the root node of the index, and thecurrent key to the first partial key of the root node. The method maythen repeat the steps of: (1) comparing the partial key of the searchvalue to the current key; (2) when the search partial key is less thanthe current key, setting the current node to a node identified based onan offset added to the child group pointer of the current node andsetting the current key to the first partial key of the identified node;(3) when the search partial key is greater than the current key, settingthe current key to a next partial key of the current node, and (4) whenthe search partial key is equal to the current key, comparing the searchvalue with a record value identified by the full key pointercorresponding to the current key. The method may involve repeating thesteps until locating the search value in a leaf node or determining thatno leaf node corresponds to the search value.

Consistent with other disclosed embodiments, a system for locating asearch value in a database. The system may comprise a cache memoryhaving a cache line size and storing an index comprising groups of nodesthat comprise a child group pointer, a number of fixed size partialkeys, and a number of full key pointers corresponding to the number offixed size partial keys. The system may also comprise a processor; and amemory storing instructions that, when executed by the processor, causethe processor to perform operations. The operations may comprisecomputing a partial key for the search value and setting a current nodeto the root node and setting a current key to the first partial key ofthe root node. The operations may further comprise repeating the stepsof: (1) comparing the partial key of the search value to the currentkey; (2) when the search partial key is less than the current key,setting the current node to a node identified based on an offset addedto the child group pointer of the current node and setting the currentkey to the first partial key of the identified node; (3) when the searchpartial key is greater than the current key, setting the current key toa next partial key of the current node, and (4) when the search partialkey is equal to the current key, comparing the search value with arecord value identified by the full key pointer corresponding to thecurrent key.

Consistent with other disclosed embodiments, computer-readable media,such as storage devices, may store program instructions that areexecutable by one or more processors to implement any of the methods,disclosed herein.

It is to be understood that both the foregoing general description andthe following detailed description are exemplary and explanatory onlyand are not restrictive of the disclosed embodiments, as claimed.

BRIEF DESCRIPTION OF THE DRAWINGS

The accompanying drawings, which are incorporated in and constitute apart of this specification, illustrate several embodiments and, togetherwith the description, serve to explain the disclosed principles. In thedrawings:

FIG. 1 is a diagram illustrating a cache-sensitive index, known in theprior art;

FIG. 2 is a diagram illustrating a fixed-size partial key index, knownin the prior art;

FIG. 3 is a diagram illustrating a computer system capable ofimplementing a cache-sensitive index using partial-keys, consistent withdisclosed embodiments;

FIG. 4 is a diagram illustrating a node of a cache-sensitive index usingfixed-size partial keys, consistent with disclosed embodiments;

FIG. 5 is a diagram illustrating a node group for a cache-sensitiveindex using fixed-size partial keys, consistent with disclosedembodiments;

FIG. 6 is a diagram illustrating a tree representation of acache-sensitive index using fixed-size partial keys, consistent withdisclosed embodiments;

FIG. 7 is a diagram illustrating an example of a cache-sensitive indexusing fixed-size partial keys populated with index data, consistent withdisclosed embodiments;

FIG. 8 is a flow diagram illustrating an exemplary process forbulk-loading a cache-sensitive index using fixed-size partial keys,consistent with disclosed embodiments;

FIG. 9 is a pseudo-code example of a process for creating a sorted listof database records for the bulk-loading process, consistent withdisclosed embodiments;

FIG. 10 is a flow diagram illustrating an exemplary process for creatingnon-leaf nodes during the bulk-load process, consistent with disclosedembodiments;

FIGS. 11A and 11B are a pseudo-code example of an ancestral key-copyingprocess for creating non-leaf nodes during the bulk-load process,consistent with disclosed embodiments; and

FIGS. 12A and 12B are flow diagrams illustrating an exemplary deletionprocess of a cache-sensitive index using fixed-size partial keys,consistent with disclosed embodiments.

DESCRIPTION OF THE EMBODIMENTS

Disclosed embodiments provide methods and systems for implementing acache-sensitive index that also uses partial-keys to improve data lookuptimes. A node in such an index include one pointer, as with acache-sensitive B+ tree. In addition, the index includes partial keys tospeed up comparison. In some embodiments, the partial keys of each nodeare configured to fit within a cache line, such as 64 bytes, so thattraversing the partial keys does not require crossing a cache line. Thisalso improves the look-up speed of the index.

In certain embodiments, the format of the partial key may be variable,such that the offset may take one or two bytes depending on the value ofthe offset. Because of the addition of partial keys to thecache-sensitive structure, prior solutions for deleting records from acache-sensitive index by marking deleted records as deleted cannot beused. Thus, a method is provided for maintaining live updates of thecache-sensitive partial-key index, providing for live, physical deletionof records, and a bulk-loading method for a full reorganization of theindex.

Reference will now be made in detail to exemplary embodiments, examplesof which are illustrated in the accompanying drawings. Whereverconvenient, the same reference numbers will be used throughout thedrawings to refer to the same or like parts.

FIG. 3 is a diagram illustrating a computer system 300 capable ofimplementing disclosed embodiments, including exemplary systemcomponents. The components and arrangement, however, may be varied.Computer system 300 may include a processor 305, a memory 310,input/output (I/O) devices 330, cache memory 325, and storage 320.Computer system 300 may be implemented in various ways. For example,computer system 300 may be a general purpose computer, a server, amainframe computer, or any combination of these components. Computersystem 300 may communicate over a link with a network (not shown). Forexample, the link may be a direct communication link, a LAN, a WAN, orother suitable connection. The network may include the Internet.Computer system 300 may be standalone or it may be part of a subsystem,which may, in turn, be part of a larger system, such as a legacy nameserver system.

Processor 305 may include one or more known processing devices, such asa microprocessor from the Pentium™ or Xeon™ family manufactured byIntel™, the Turion™ family manufactured by AMD™, or any of variousprocessors manufactured by Sun Microsystems. Memory 310 may include oneor more storage devices, including main-memory devices, configured tostore information used by processor 305 to perform certain functionsrelated to disclosed embodiments. Storage 320 may include a volatile ornon-volatile, magnetic, semiconductor, tape, optical, removable,nonremovable, or other type of storage device or computer-readablemedium, including mass storage devices. Cache memory 325 may include L1cache on the same chip as processor 302, static RAM, or dynamic RAM.Computer system 300 may store databases, as memory space permits, inwhole or in part in a main memory system such as memory 310.

In one embodiment, memory 310 may include one or more index maintenanceprograms or subprograms 315 loaded from storage 320 or elsewhere that,when executed by central repository server 120, perform variousprocedures, operations, or processes consistent with disclosedembodiments. For example, memory 310 may include a bulk loading programthat periodically rebuilds the entire cache-sensitive, partial-keyindex; a deletion program that facilitates live deletions from thecache-sensitive, partial-key index; an insertion program thatfacilitates live additions to the cache-sensitive, partial-key index;and an integrative support program that links the other programs andallows them to use a common database, provides a common user interfacefor setting system parameters, performs basic bookkeeping tasks, andprovides guidance and help. Memory 310 may also include other programsthat perform other functions and processes, such as programs thatprovide communication support, Internet access, etc.

Methods, systems, and articles of manufacture consistent with disclosedembodiments are not limited to separate programs or computers configuredto perform dedicated tasks. For example, memory 310 may be configuredwith a index maintenance program 315 that performs several functionswhen executed by processor 305. For example, memory 310 may include asingle program 315 that performs the index maintenance functions, orprogram 315 could comprise multiple programs. Moreover, processor 305may execute one or more programs located remotely from computer system300. For example, computer system 300 may access one or more remoteprograms that, when executed, perform functions related to disclosedembodiments.

Memory 310 may be also be configured with an operating system (notshown) that performs several functions well known in the art whenexecuted by processor 305. By way of example, the operating system maybe Microsoft Windows™, Unix™, Linux™, Solaris™, or some other operatingsystem. The choice of operating system, and even to the use of anoperating system, is not critical to any embodiment.

Computer system 300 may include one or more I/O devices 330 that allowdata to be received and/or transmitted by computer system 300. I/O 330devices may also include one or more digital and/or analog communicationinput/output devices that allow computer system 300 to communicate withother machines and devices, such as a user's computing device or someother a client computer (not shown). Client computers may providerequests from users representing queries of the data stored in storage320 or memory 310. Computer system 300 may receive data from externalmachines and devices and output data to external machines and devicesvia I/O devices 330. The configuration and number of input and/or outputdevices incorporated in I/O devices 330 may vary as appropriate forcertain embodiments.

Computer system 300 may include one or more databases that storeinformation and are access and/or managed through computer system 300.The databases may be stored in storage 320 or memory 310. The databasesmay include, for example, data and information related to domain names,IP addresses, and other information needed to process a DNSSEC request.In some embodiments, the databases or other files may include datasimilar to the items shown in FIG. 4. Systems and methods of disclosedembodiments, however, are not limited to separate databases.

FIG. 4 is a diagram illustrating a node 400 for a cache-sensitive indexusing fixed-size partial keys, consistent with disclosed embodiments.Node 400 may be a root, intermediate, or leaf node. Node 400 may includechild group pointer 405. Child group pointer 405 may be a downlinkpointer to a child group. As discussed above, a group is a collection ofnodes stored contiguously in memory. Child group pointer 405 may containthe address of the first node in the group of nodes. The location of theremaining nodes may be calculated by adding the size of each node to theaddress of the pointer.

Node 400 may also contain partial keys 410. In certain embodiments,partial keys 410 comprise 4 bytes, the 4 bytes representing two distinctpieces of information: 1) the starting position from the base key wherethe bytes begin to differ (i.e. the offset); and 2) the first two orthree bytes of the key itself (key data). Although fixed in size, thenumber of bytes used by the offset and the key data may vary. Forexample, if the starting position from the base key occurs within thefirst 254 bytes, then the offset may use one byte and the key data mayuse three bytes of the partial key. On the other hand, when the startingposition from the base key occurs at or beyond the 254^(th) position,the offset may use two bytes and the key data may use two bytes. Thevariable format of the partial key allows the key data to consume morebytes in most cases, decreasing the amount of time a search key willmatch exactly a partial key. Traversing the index using the partial keysis explained in more detail with regard to FIG. 7 below.

In some embodiments all partial keys 410 of node 400 may be stored inthe first block of node 400. In other words, child group pointer 405 andpartial keys 410 fit within one cache line. This allows faster traversalof the partial keys because there is no possibility of a cache miss.Although FIG. 4 shows 15 partial keys 410 per node 400, any number ofpartial keys may be used, as long as the number of partial keys 410 fitin the first cache line. Therefore, the larger the cache line, the morepartial keys 410 that may fit in a node 400. A smaller cache line mayrequire that node 400 contain fewer partial keys 410 than shown in FIG.4.

Node 400 may also include full key pointers 415. Full key pointers maybe stored as part of node 400 for situations where the partial keycompletely matches the data that is the subject of a search. Full keypointers 415 point to a record (row) in the database, and make the fullkey available for comparison should a partial key comparison fail toresolve the query. The number of full key pointers 415 will match thenumber of partial keys 410 per node. Node 400 may also include recordcount 420. Record count 420 may indicate the number of partial keys 410and full key pointers 415 that have data (i.e. are not null) andprocessor 305 may use record count 420 during a record insertion,deletion, or bulk load process.

The size of child group pointer 405 and full key pointers 415 may varydepending on the size of the database. For example, child group pointer405 and full key pointers 415 may typically be four bytes long. In otherembodiments, child group pointer 405 and full key pointers 415 may beeight bytes long, allowing for a larger maximum size for the database.The size of the pointers is not important, so long as enough bytes areused to accommodate the maximum size of the database and the size isconsistent between all child group pointers 405 and full key pointers415 within the index.

FIG. 5 is a diagram illustrating a node group 500, consistent withdisclosed embodiments. Node group 500 is a collection of nodes 400organized contiguously in memory to allow sub-tree pointers to berepresented by child group pointer 405. Node group 500 may also includea node count that indicates the number of nodes 400 in the group thatare populated with data (i.e. are not null). Processor 305 may use thenode count during record insertion and deletion processes or the bulkloading process.

FIG. 6 is a diagram illustrating a tree representation of acache-sensitive index 600 using partial keys, consistent with disclosedembodiments. Index 600 of FIG. 6 may represent an index tree for adatabase that contains 1,024 records. Each record in the database may berepresented by a pointer in one of the leaf nodes. In the example ofFIG. 6, Groups 6-68 contain the leaf nodes and Groups 1-5 containintermediate nodes. Child group pointers 405 of each node in Groups 2-5may contain the address of a group of leaf nodes. Child group pointer405 of each node in Group 1 may contain the address of a group ofintermediate nodes, e.g. Groups 2-5. The root of the tree may containone node. The child group pointer 405 of the root node may point to onegroup, e.g. Group 1.

The number of nodes 400 that have data in Group 1 correspond to thenumber of partial keys 410 and full key pointers 415 that have data inthe root node. In some embodiments, the number of nodes in the childgroup may be one higher than the number of keys in a node. In suchembodiments, the number of nodes populated in the child group will beone higher than the number of keys populated. For example, if all keysin the root node are populated with data, then the tree may have sixteennodes 400 in Group 1 populated with data. In such embodiments if asearch key is greater than the last key in a root or intermediate node,then the next node searched is the rightmost node in the child group. Inother embodiments, the number of nodes in the child group may equal thenumber of keys in a node. In the example of FIG. 6, which has one morenode in each group than keys in the node, the root node has three keyspopulated so it's child group, Group 1, has four nodes populated.Similarly, Node 4 of Group 1 has 14 keys populated with data so itschild group, Group 5, has 15 nodes 400 populated with data. Node 1 hasall 15 keys populated with data so its child group, Group 2, has 16nodes populated with data.

An example of traversing index 600 will now be explained using FIG. 7.Index 600 may be stored, for example, in cache memory 325 of system 300and traversed using processor 305. In the example of FIG. 7, index 600contains five keys per node and uses 4 byte addressing, so each noderequires 48 bytes. Processor 305 may receive a search key andinstructions to determine whether a record exists in a database thatmatches the search key. For example, processor 305 may receive thesearch key “Abner.” Processor 305 may compute a partial search key, forexample “0Abn,” from the search key. For readability, FIG. 7 shows thepartial key values in ASCII. However, to fit the three ASCII charactersinto four bytes, the data would be encoded as hexadecimal values. Forexample, “0Abn” may actually be stored as 0x0041626E.

Processor 305 may then compare the search key with the first partial key410 of the root node of the index. In the example of FIG. 7, the Node 1is the root node, and the first partial key contains “0Abb.” Thus,processor 305 may compare the three letters of the search key startingat position zero with partial key 1.

Processor 305 may determine that the search key is greater than partialkey 1 in the root node and, therefore, perform a comparison with partialkey 2. When a search key is greater than a partial key, processor 305may determine that a comparison with the next partial key in the node isrequired. To recompute the partial search key for comparison with thenext partial key in the node, processor 305 may determine whether thefirst three bytes of the partial search key match the first three bytesof partial key 1 of the root node. If so, processor 305 may increase theoffset of the partial search key by 2. If the first three bytes do notmatch, processor 305 may determine if the first two bytes match. If so,processor 305 may increase the offset by one. If only the first bytematches, which represents the offset, then processor 305 may not make anadjustment to the search partial key. In embodiments with a variablesize partial key, when the increase to the offset causes the offset toexceed a predetermined limit, such as 250, processor 305 may adjust thekey data portion of the partial key to use two bytes instead of three.Processor 305 may then adjust the key data in the partial search keyaccording to the new offset, if needed. In the example of FIG. 7,processor 305 may determine that the first three bytes match (the threebytes including the offset), and add two to the offset, resulting in anew partial search key of “2ner.”

Processor 305 may compare the partial search key with partial key 2 ofthe root node. Processor 305 may determine that the search key is lessthan partial key 2. When a search key is less than a partial key, theprocessor may determine that the next node to search corresponds withthe child node of the partial key. For example, the child node ofpartial key 2 of the root node is the second node of the groupidentified by the child group pointer 405 of the root node. Processor305 may determine the address of the child node by adding an offset tothe value of child group pointer in Node 1. For example, because eachnode is 48 bytes long, the offset may be a multiple of 48.

Processor 305 may determine the amount of the offset by multiplying theabsolute position of the partial key within the node by the node size.For example, the absolute position of the first partial key in a node iszero, the absolute position of the second partial key is one, etc. Inthe present example, processor 305 may determine that the offset is 48because the absolute position of partial key 2 of the root node is one.Thus, processor 305 may calculate the address of the second node in thechild group by adding 48 to the value of the child group pointer.Similarly, if processor 305 used partial key 4 of the root node in thecomparison, processor 305 may determine that the offset is 144 (i.e. theabsolute position of three multiplied by the node size of 48).

After calculating the address of the second node of the child group, andsetting the current node to the second node of the child group,processor 305 may compare the value of the partial search key with thevalue of the first partial key of the current node, e.g. comparing“2ner” with “2iga.” When processor 305 determines that the search key isgreater than the first partial key, processor 305 may recompute thepartial search key as discussed above. In the present example, processor305 may not make any adjustment because only the first byte of thepartial search key matches partial key 1 of the current node.

Processor 305 may then compare the partial search key with the secondpartial key of the current node, e.g. comparing “2ner” with “21e.” Whenprocessor 305 determines that the search key is greater than the secondpartial key, processor 305 may determine that no adjustment is necessaryto the partial key (because only the offset matches) and proceed tocompare the partial search key with the third partial key of the currentnode. Because the partial keys match, processor 305 may perform a fullkey comparison, by retrieving the value of the full key. Processor 305may retrieve the value “Abner,” which is stored at the addressrepresented by the third full key pointer of the current node. Becausethe search key matches the retrieved data, Processor 305 may determinethat a record has been found, and may return the location of theretrieved record.

This example demonstrates how processor 305 never needed to do anyexpensive full key comparisons until the leaf node was found. All of theinformation required to traverse the index and determine the next childnode is contained in partial keys 410.

In another example, processor 305 may receive a search key of “Abramo.”As described above, processor 305 may compute a partial search key andcompare the partial search key with the first partial key of the rootnode, e.g., comparing “0Abr” with “0Abb.” Because the value of thesearch key is greater than the value of the first partial key, processor305 may recompute the partial search key, adding two to the offset andadjusting the key data. Processor 305 may then compare the partialsearch key “2ram” to the second partial key of the root node “2ram.”Because the search key matches the partial key, processor 305 mustperform a full key comparison by retrieving the value of the dataresiding at the address in the second full key pointer. Processor 305retrieves the record, e.g. “Abram,” and compares “Abram” with “Abramo.”Because the full key comparison shows the search key is greater,processor would continue with the next partial key. In the example ofFIG. 7, the next partial key is null (empty).

In disclosed embodiments where the number of nodes in a group is onemore than the number of keys in a node, this indicates that if Abramoexists in the tree, it will be in the right subtree (represented in FIG.7 by the ellipsis after the second node). Processor 305 may not need tostore a partial key or full key pointer in the root or intermediatenodes for the right subtree, as all searches beyond the last populatedkey must traverse down the right subtree. In other embodiments where thenumber of nodes in a group equals the number of keys in a node, if thenext partial key is null (empty), this results in a determination thatthe search key does not exist in the database. Regardless of theembodiment used, if the next partial key did contain data, processor 305would continue its search until either a record is found, or theprocessor determines that no leaf node matches the search key (i.e. norecord exists in the database that matches the search key).

As discussed above, partial keys 410 may be encoded in a total byteordered representation. This puts the most significant bytes in thebeginning of the key and increases the speed of comparisons. However,data in the rows of the database may not be stored that way. In such anembodiment, processor 305 may need to dereference the full key before acomparison of the search key and the full key.

FIG. 8 is a flow diagram illustrating an exemplary load process 800 fora cache-sensitive index using fixed size partial keys. Processor 305 mayperform process 800 to rebuild the index as part of an operationalprocedure. A bulk load builds the index with completely full nodes toprovide optimal read performance. For example, a B+ tree performs betterwith the nodes are full and balanced. In certain embodiments, process800 may be implemented according to index maintenance programs 315 inmemory 310. Process 800 may build all of the leaf nodes in a totalsorted order first before building the higher level nodes in the tree.The leaf nodes may be sorted using a non-recursive merge sort that usesthe partial keys, decreasing the load time.

In step 805, processor 305 reads a record from an input, such as a file,and populates the first key in the first node of a group. For example,as discussed with respect to FIGS. 4 and 5, in the first node of thegroup, the child group pointer 405 of the group is null (empty), and thefirst full key pointer 415 points to the record that processor 305 justread. The first partial key may contain an offset of zero and the firstthree bytes of the key of the record. Processor 305 may encode thepartial key in a total byte order representation (big Endean). In step810, processor 305 may push this group onto a stack. The stack storesgroups, with the most recently added group at the top of the stack.Further, each group contains nodes, and each node contains a number ofslots (partial keys and full key pointers). Each full key pointerrepresents one record.

In step 815, processor 305 determines whether the number of records(i.e. populated keys in the nodes of the group) in the top two stackentries form a power of two (e.g., 2, 4, 6, 8, 16, 32, . . . ). If thenumber of records does form a power of 2 (step 815, Yes), then processor305 may pop the two stack entries and combine them into a new sortedlist. For example, if the top two stack entries contain a group withthree records (e.g., the first three keys of the first node arepopulated), and a group with one record (e.g., the first key of thefirst node is populated), then processor 305 may combine the two groupsinto a single, sorted group and push the newly sorted group onto thestack. FIG. 9 depicts pseudo-code for combining two entries into a newsorted list that may be used in some embodiments.

In some situations, the number of entries may fit in one group. In thissituation, processor 305 may create one group for the sorted list.However, if the number of entries do not fit in one group, processor 305may create the sorted list by linking groups of nodes. For example, ifprocessor 305 combines a group with sixteen populated nodes (i.e. 240records, if each node stores 15 keys and each group contains 16 nodes)and a group with three populated nodes, the sorted list that resultsfrom the combination may comprise two groups of nodes, with the childgroup pointer of the first group containing the address of the secondgroup. After this manner, the sorted list may contain a linked list ofgroups of nodes.

After processor 305 creates the new sorted list then, in step 825, thenew sorted list is pushed onto the stack. The process then continues atstep 815, where processor 305 checks to see if the number of records inthe new sorted list and the number of records in the stack entry belowthe new sorted list form a power of two. If they do, processor 305repeats steps 820 and 825. If the top two stack entries do not form apower of two then, in step 830, processor 305 determines if there areany more records left to read in the database. If so (step 830, No),processor 305 repeats steps 805-825 as needed.

If all records have been read (step 830, Yes) then, in step 835,processor 305 may combine all the groups on the stack, as discussed withregard to step 820, resulting in one sorted, linked-list of groups.Processor 305 may also balance the two groups at the end of the sortedlist. For example, the final node may be less than half full of keys andthe final group may be less than half full of nodes. When this occurs,processor 305 may rebalance the final nodes of the final group, or thefinal two groups of the linked list so that each group is at least halffull of nodes, and each node is at least half full of keys. Aftercompleting step 835, processor 305 has created groups of nodes that mayserve as leaf nodes in the cache-sensitive index using fixed sizepartial keys. In step 840, processor 305 creates the non-leaf nodes,such as the root and intermediate nodes.

FIG. 10 is a flow diagram of a process 1000 that creates the non-leafnodes during a bulk-loading process, consistent with disclosedembodiments. Processor 305 may perform process 1000 as part of step 840of FIG. 8. In certain embodiments, process 1000 may be implementedaccording to index maintenance programs 315 in memory 310.

In step 1005, processor 305 may calculate the total nodes by countingthe number of populated nodes in the sorted list of leaf groups.Processor 305 may have created the sorted list of leaf groups as part ofprocess 800. In step 1010, processor 305 may calculate the number ofgroups in the next higher level by dividing the total nodes by thenumber of nodes in a group, rounding to the next highest number. Forexample, if the sorted list of leaf nodes contains 63 nodes, and eachgroup contains 16 nodes, then processor 305 may determine that the levelabove the leaf nodes requires four groups (63/16=3.9, rounded up to 4).However, if processor 305 counts 1983 leaf nodes, then processor 305 maycreate 124 groups. In step 1015, processor 305 may create the groups inthe next higher level by allocating memory for the groups and creating alinked list of the groups by setting the child group pointer of thefirst group to the address of the second group etc. Processor 305 mayset all other values and pointers in the groups to zero (null). Forpurposes of discussion, this level may be L1 (the level above the leafnodes).

Next, in step 1020, processor 305 determines whether the calculatednumber of groups (i.e. the number of groups just created) is less thanthe number of nodes in a group. If not (step 1020, No), processor 305repeats steps 1010 and 1015 for higher layers in the tree. For example,using the example of 1983 leaf nodes discussed above, processor 305created 124 groups in L1. Because 124 is greater than 16, processor 305may repeat steps 1010-1020 for a higher layer, for example L2 (the levelabove L1). In this example, processor 305 may create 8 group nodes inL2, because 124/16=7.75, which is rounded up to 8.

If the calculated number of groups is less than the number of nodes in agroup (step 1020, Yes) then, in step 1025, processor 305 creates a rootgroup with one node. At this point, processor 305 has created allrequired groups for the index and must now populated the keys, partialkeys, and down pointers of the non-leaf nodes. To accomplish this, instep 1030, processor 305 may iterate all of the leaf nodes and copy thekeys and partial keys from the rightmost key in the leaf node up to thenext highest non-leaf node. As processor 305 populates the first key ineach node, processor 305 may set the child group pointer to the addressof the group below that supplied the key value. FIG. 11A depictspseudo-code for the leaf-node iteration that may be used in someembodiments.

In step 1035, processor 305 populates the nodes and groups of the higherlevel intermediate nodes, again copying the rightmost key in each nodeto next highest node, setting the child group pointers as describedabove with regard to step 1030. Processor 305 may repeat step 1035 untilthe keys of all intermediate levels, up to and including the root, havebeen populated. FIG. 11B depicts pseudo-code for the non-leaf nodeiteration that may be used in some embodiments to perform the process ofstep 1035.

In step 1040, processor 305 may set all the child group pointers in theleaf groups to null, thereby breaking the links between the leaf nodes.When processor 305 traverses the index, a child group pointer of nullindicates that the node is a leaf node and no further traversal ispossible. Therefore, if processor 305 has not found a record in theindex matching the search key and the child group pointer of the node isnull, then no matching record exists.

FIGS. 12A and 12B are flow diagrams illustrating an exemplary deletionprocess 1200 of a cache-sensitive index using fixed-size partial keys,consistent with disclosed embodiments. One challenge of maintaining apartial-key B+ tree is removal of deleted entries from the index. Priorsolutions use a lazy deletion process that flags deleted entries in theindex, leaving the node structure untouched until a clean-up process,such as a bulk load, is performed. Lazy deletion, however, leads tolookup speed degradation and poor space usage in a cache-sensitiveindex. Process 1200 offers a live-update solution for the deletionproblem. Before process 1200 begins, processor 305 may receive a key tobe deleted. This may be known as the target key. In certain embodiments,process 1200 may be implemented according to index maintenance programs315 in memory 310.

At Step 1205, processor 305 finds the group (i.e. the target group) andthe leaf node (i.e. the target node) that store the key to be deleted.Processor 305 may find the target node by traversing the index asdiscussed above with regard to FIG. 7. Once processor 305 finds thetarget node, in step 1210 processor 305 determines whether the targetnodes has more than the minimum number of keys. In a B+ tree structure,each node except the root node should be at least half full. Thus, inthe example of a node having 15 keys, the minimum number of keys isseven. If the target node does not have the minimum number of keys (step1210, No), then processing continues at 1255, which will be discussed inmore detail with regard to FIG. 12B.

If the target node does have more than the minimum number of keys (step1210, Yes), then in step 1215 processor 305 determines whether thetarget key occupies the last slot of the target node. As used herein, aslot represents a partial key and its corresponding full key pointer. Ifthe target key is not in the last slot (step 1215, No) then, in step1220 processor 305 may remove the target key from the target node andshift the keys to the right of the target key one slot to the left.After shifting the keys to the left, processor 305 may re-compute thepartial key of the slot that had been immediately to the right of thetarget key. Processor 305 has then completed the deletion, and process1200 ends.

If the target key occupies the last slot in the target node (step 1215,Yes) then, in step 1225, processor 305 may remove the target key fromthe target node. This causes the formerly penultimate slot to become thenew last key for the target node. After removing the target key, in step1230, processor 305 may determine whether the target node is the lastnode of the target group. The last node in a group is the rightmost nodethat contains data. For example, if a group can have 15 nodes, but fiveof the nodes are empty, then the tenth node is the last node of thegroup.

If the node is not the last node of the target group (step 1230, No),then in step 1235 processor 305 updates the keys in the parent node andupdates the partial key of the first slot in the node to the right ofthe target node (i.e. the right sibling node). Processor 305 may updatethe partial key of the right sibling node to reflect a new offset andnew key data when compared with the new last key (in the formerlypenultimate slot). Furthermore, because the slots in the parent nodecontain the value of the key in the last slot of each child node,processor 305 must copy the new last key of the target node to thecorresponding slot in the parent node, adjusting the partial key of theupdated slot in the parent and the partial key to the right of theupdated slot in the parent node. Processor 305 has then completed thedeletion and process 1200 ends.

If the node is the last node of the target group (step 1230, Yes) then,in step 1240, processor 305 updates the partial key of the first slot ofthe group to the right of the target group (i.e. the right adjacentgroup or “RAG”). If there is no right adjacent group (i.e. the deletedkey is the highest key of the tree), then processor 305 has finished thedelete and process 1200 ends. The partial key of the RAG will beadjusted in the same manner as described with regard to the rightsibling node, as described above with respect to step 1235.

In step 1245, processor 305 may locate the common node ancestor (“CNA”)of the target node and the RAG. The common node ancestor is the lastnode in the path from the root node to the leaf node that is common toboth the target node and the first node of the RAG. Once processor 305has located the CNA, processor 305 may update the slot in the CNA thatcontains the target key (because the target key was the last slot in thetarget node). Processor 305 may use the formerly penultimate slot of thetarget to update the slot in the CNA. After updating the slot in theCNA, processor 305 may then re-compute the partial key of the slot inthe CNA immediately to the right of the updated slot. In step 1250,processor 305 may update the partial keys of all the nodes traversedwhen traversing the tree from the updated slot of the CNA to the RAG.Processor 305 has then completed the deletion and process 1200 ends.

FIG. 12B represents process 1255, which processor 305 performs when thetarget node does not have more than the minimum number of keys (step1210, No). To ensure the index tree is balanced, the target node cannotbe left with less than the minimum number of keys after deletion of thetarget key. Because of this, the nodes must be rebalanced after deletionof the target key. In step 1260, processor 305 determines if any node inthe target group contains more than the minimum number of slots. If atleast one other node in the target group has more than the minimumnumber of slots (step 1260, Yes) then, in step 1265, processor 305redistributes the slots evenly across all nodes of the target group andupdates the parent node according to the new key distribution. Processor305 may also need to update grandparent nodes to reflect the newdistribution. If the target key was the last key of the target group(step 1267, Yes), then processor 305 will perform steps 1240 to 1250, asdescribed above with regard to FIG. 12A, to adjust the keys in the RAGand CNA. If the target key was not the last key of the target group(step 1267, No), processor 305 has completed the deletion and process1200 ends.

If no nodes in the target group contain more than the minimum number ofkeys then, in step 1270, processor 305 migrates the target node's slotsto an adjacent node of the target group and deletes the target node.Because none of the group's nodes contain more than the minimum number,any node in the group can fit the slots of the target node. Once thetarget node is deleted, processor 305 determines whether the targetgroup now contains less nodes than the minimum required nodes. Forexample, when a node may have a maximum of 15 keys and a minimum of 7keys, each group may have at most 16 nodes and at least 8 nodes. If thenumber of nodes falls below the minimum number, the nodes of theadjacent groups should be redistributed to keep the tree balanced.

If the target group, after deletion of the target node, contains morenodes than the minimum required (step 1275, No) then, processor 305 hasfinished the deletion and process 1200 ends. If the target groupcontains less than the minimum number of nodes after deletion of thetarget node (step 1275, Yes) then processor 305 determines whether anadjacent sibling group (i.e. “sibling groups” immediately to the left orto the right) contains more than the minimum number of nodes. If one ofthe sibling groups does contain more than the minimum number of nodes(step 1280, Yes) then, at step 1285, processor 305 moves one of thenodes from the sibling group into the target group. Processor 305 thenupdates the two slots in the parent node and potentially slots in thegrandparent node. After updating all appropriate slots then, in step1287, processor 305 determines whether the target key was the rightmostkey of the target group and if so, whether the migrated node came fromthe left sibling group. If both answers are yes (step 1287, Yes), thenprocessor 305 performs steps 1240-1250, as explained above. If not (step1287, No), then processor 305 has finished the deletion and process 1200ends.

If none of the sibling groups contain more than the minimum number ofnodes (step 1280, No) then, in step 1290, processor merges the targetgroup with a sibling adjacent group. This causes processor 305 to deleteone node in the parent group and update the grandparent node. Aftermerging the two groups, processor 305 may determine whether the parentgroup, with the node deleted, now contains less than the minimum numberof nodes. If so (step 1295, Yes), then processor 305 may perform steps1280, 1285, 1287, 1290, and 1295, as appropriate, using the parent groupas the “target group.”

If the parent group does have the minimum number of nodes required (step1295, No), then processor 305 returns to step 1287 to determine whetherthe target group was merged with the left sibling group and the targetkey the rightmost key of the target group before deletion. If so (step1287, Yes), processor 305 repeats steps 1240-1250. If not (step 1287,No), processor 305 has finished the deletion and process 1200 ends.

While process 1200 was described as a process of removing keys fromnodes and adjusting nodes up the tree, in order to obtain transactionalbehavior, processor 305 must perform any changes to a node in a singleatomic step. Changing two items in a node, such as a partial key and afull key pointer, is non-atomic. To accommodate such non-atomic changes,processor 305 may make a copy of the leaf group, execute process 1200 tomake the changes in the copy, and swap in the new group after completingthe changes by updating the child group pointer of the parent node. Inmore complex updates, the processor may copy subtrees of the index upuntil a common ancestor and swap in the entire subtree after all changeshave been made.

The foregoing descriptions have been presented for purposes ofillustration and description. They are not exhaustive and do not limitthe disclosed embodiments to the precise form disclosed. Modificationsand variations are possible in light of the above teachings or may beacquired from practicing the disclosed embodiments. For example, thedescribed implementation includes software, but the disclosedembodiments may be implemented as a combination of hardware and softwareor in firmware. Examples of hardware include computing or processingsystems, including personal computers, servers, laptops, mainframes,micro-processors, and the like. Additionally, although disclosed aspectsare described as being stored in a memory on a computer, one skilled inthe art will appreciate that these aspects can also be stored on othertypes of computer-readable storage media, such as secondary storagedevices, like hard disks, floppy disks, a CD-ROM, USB media, DVD, orother forms of RAM or ROM.

Computer programs based on the written description and disclosed methodsare within the skill of an experienced developer. The various programsor program modules can be created using any of the techniques known toone skilled in the art or can be designed in connection with existingsoftware. For example, program sections or program modules can bedesigned in or by means of .Net Framework, .Net Compact Framework (andrelated languages, such as Visual Basic, C, etc.), XML, Java, C++,JavaScript, HTML, HTML/AJAX, Flex, Silverlight, or any other now knownor later created programming language. One or more of such softwaresections or modules can be integrated into a computer system or existingbrowser software.

Other embodiments will be apparent to those skilled in the art fromconsideration of the specification and practice of the embodimentsdisclosed herein. The recitations in the claims are to be interpretedbroadly based on the language employed in the claims and not limited toexamples described in the present specification or during theprosecution of the application, which examples are to be construednon-exclusive. Further, the steps of the disclosed methods may bemodified in any manner, including by reordering steps and/or insertingor deleting steps. It is intended, therefore, that the specification andexamples be considered as exemplary only, with a true scope and spiritbeing indicated by the following claims and their full scopeequivalents.

What is claimed is:
 1. A computer-implemented method of locating asearch value in a database, the database having an index comprisinggroups of nodes, wherein each node comprises: a child group pointer, anumber of fixed size partial keys, and a number of full key pointerscorresponding to the number of fixed size partial keys, the methodcomprising: computing a search value partial key for the search value,wherein the search value partial key is a fixed size and comprises afirst byte offset and a subset of the search value; setting a currentnode to a root node and setting a current partial key to a first partialkey of the root node, wherein the first partial key of the root node isthe fixed size and comprises the first byte offset and a subset of a keyreferenced by a first full key pointer of the root node; and comparingthe search value partial key to the current partial key, wherein thecomparing comprises; when the search value partial key is less than thecurrent partial key, setting the current node to a node referenced by achild group pointer of the current node and setting the current partialkey to a first partial key of the identified node, wherein the firstpartial key of the identified node comprises a second byte offset and asubset of a key referenced by a first full key pointer of the identifiednode, when the search value partial key is greater than the currentpartial key, setting the current partial key to a next partial key ofthe current node, wherein the next partial key of the current node isthe fixed size and comprises a third byte offset and a subset of a keyreferenced by a next full key pointer of the current node, when thesearch value partial key is equal to the current partial key, comparingthe search value with a key referenced by a full key pointercorresponding to the current partial key, and determining that thesearch value matches the key referenced by the full key pointer; andoutputting a record value corresponding to the key referenced by thefull key pointer in response to the determining.
 2. Thecomputer-implemented method of claim 1, wherein the node referenced bythe child group pointer is identified based on an offset and the offsetcomprises a node size multiplied by a position of the current partialkey within the current node.
 3. The computer-implemented method of claim1, wherein when the search value is greater than the key referenced bythe full key pointer corresponding to the current partial key, themethod further comprises performing the setting of the current partialkey to the next partial key of the current node.
 4. Thecomputer-implemented method of claim 1, wherein when the search value isless than the key referenced by the full key pointer corresponding tothe current partial key, the method further comprises performing thesetting of the current node to the node referenced by the child grouppointer of the current node and the setting of the current partial keyto the first partial key of the identified node.
 5. Thecomputer-implemented method of claim 1, wherein when the search valuepartial key is greater than the current partial key, the method furthercomprises recalculating the search value partial key.
 6. Thecomputer-implemented method of claim 5, wherein recalculating the searchvalue partial key comprises: calculating a new byte offset by: addingtwo to the first byte offset when the first three bytes of the searchvalue partial key match the first three bytes of the current partialkey, and adding one to the first byte offset when the first two bytes ofthe search value partial key match the first two bytes of the currentpartial key; and adjusting the subset of the search value according tothe new byte offset.
 7. The computer-implemented method of claim 6,wherein when the new byte offset exceeds a predetermined value, the newbyte offset comprises an additional byte and the subset of the searchvalue comprises one less byte.
 8. The computer-implemented method ofclaim 1 further comprising: determining that the search value is notlocated in the database based on the comparing the search value partialkey to the current partial key; and outputting an indication that thesearch value is not located in the database in response to thedetermining.
 9. A computer readable medium storing instructions forcausing a processor to perform the computer-implemented method ofclaim
 1. 10. A system for locating a search value in a database,comprising: a cache memory having a cache line size and storing an indexcomprising groups of nodes, wherein each node of the groups of nodescomprises: a child group pointer, a number of fixed size partial keys,and a number of full key pointers corresponding to the number of fixedsize partial keys; a processor; and a memory storing instructions that,when executed by the processor, cause the processor to performoperations comprising: computing a search value partial key for thesearch value, wherein the search value partial key is a fixed size andcomprises a first byte offset and a subset of the search value; settinga current node to a root node and setting a current partial key to afirst partial key of the root node, wherein the first partial key of theroot node is the fixed size and comprises the first byte offset and asubset of a key referenced by a first full key pointer of the root node;and comparing the search value partial key to the current partial key,wherein the comparing comprises; when the search value partial key isless than the current partial key, setting the current node to a nodereferenced by a child group pointer of the current node and setting thecurrent partial key to a first partial key of the identified node,wherein the first partial key of the identified node comprises a secondbyte offset and a subset of a key referenced by a first full key pointerof the identified node, when the search value partial key is greaterthan the current partial key, setting the current partial key to a nextpartial key of the current node, wherein the next partial key of thecurrent node is the fixed size and comprises a third byte offset and asubset of a key referenced by a next full key pointer of the currentnode, when the search value partial key is equal to the current partialkey, comparing the search value with a key referenced by a full keypointer corresponding to the current partial key, and determining thatthe search value matches the key referenced by the full key pointer; andoutputting a record value corresponding to the key referenced by thefull key pointer in response to the determining.
 11. The system of claim10, wherein the child group pointer and the number of fixed size partialkeys fit in a cache line of the cache line size.
 12. The system of claim10, wherein nodes of each group of nodes are stored contiguously in thecache memory.
 13. The system of claim 12, wherein the node referenced bythe child group pointer is identified based on an offset and the offsetcomprises a node size multiplied by a position of the current partialkey within the current node.
 14. The system of claim 10, wherein whenthe search value is greater than the key referenced by the full keypointer corresponding to the current partial key, the operations furthercomprise performing the setting of the current partial key to the nextpartial key of the current node.
 15. The system of claim 10, whereinwhen the search value is less than the key referenced by the full keypointer corresponding to the current partial key, the operations furthercomprise performing the setting of the current node to the nodereferenced by the child group pointer of the current node and thesetting of the current partial key to the first partial key of theidentified node.
 16. The system of claim 10, wherein when the searchvalue partial key is greater than the current partial key, theoperations further comprise recalculating the search value partial key.17. The system of claim 16, wherein the operation of recalculating thesearch value partial key comprises: calculating a new byte offset by:adding two to the first byte offset when the first three bytes of thesearch value partial key match the first three bytes of the currentpartial key, and adding one to the first byte offset when the first twobytes of the search value partial key match the first two bytes of thecurrent partial key; and adjusting the subset of the search valueaccording to the new byte offset.
 18. The system of claim 17, whereinwhen the new byte offset exceeds a predetermined value, the new byteoffset comprises an additional byte and the subset of the search valuecomprises one less byte.
 19. The system of claim 10, wherein each groupof nodes of the groups of nodes comprises a number of nodes, the numberof nodes being one more than the number of fixed size partial keys. 20.The system of claim 10, the operations further comprising: determiningthat the search value is not located in the database based on thecomparing the search value partial key to the current partial key; andoutputting an indication that the search value is not located in thedatabase in response to the determining.
 21. A non-transitorycomputer-readable storage medium for locating a search value in adatabase, comprising: a cache memory having a cache line size andstoring an index comprising groups of nodes, wherein each node of thegroups of nodes comprises: a child group pointer, a number of fixed sizepartial keys, and a number of full key pointers corresponding to thenumber of fixed size partial keys; and instructions which, when executedon a processor, perform a method comprising: computing a search valuepartial key for the search value, wherein the search value partial keyis a fixed size and comprises a first byte offset and a subset of thesearch value; setting a current node to a root node and setting acurrent partial key to a first partial key of the root node, wherein thefirst partial key of the root node is the fixed size and comprises thefirst byte offset and a subset of a key referenced by a first full keypointer of the root node; and comparing the search value partial key tothe current partial key, wherein the comparing comprises; when thesearch value partial key is less than the current partial key, settingthe current node to a node referenced by a child group pointer of thecurrent node and setting the current partial key to a first partial keyof the identified node, wherein the first partial key of the identifiednode comprises a second byte offset and a subset of a key referenced bya first full key pointer of the identified node, when the search valuepartial key is greater than the current partial key, setting the currentpartial key to a next partial key of the current node, wherein the nextpartial key of the current node is the fixed size and comprises a thirdbyte offset and a subset of a key referenced by a next full key pointerof the current node, when the search value partial key is equal to thecurrent partial key, comparing the search value with a key referenced bya full key pointer corresponding to the current partial key, anddetermining that the search value matches the key referenced by the fullkey pointer; and outputting a record value corresponding to the keyreferenced by the full key pointer in response to the determining. 22.The non-transitory computer-readable storage medium of claim 21, whereinthe node referenced by the child group pointer is identified based on anoffset and the offset comprises a node size multiplied by a position ofthe current partial key within the current node.
 23. The non-transitorycomputer-readable storage medium of claim 21, the method furthercomprising: determining that the search value is not located in thedatabase based on the comparing the search value partial key to thecurrent partial key; and outputting an indication that the search valueis not located in the database in response to the determining.
 24. Acomputer-implemented method of locating a search value in a database,the database having an index comprising groups of nodes, wherein eachnode comprises: a child group pointer, a number of fixed size partialkeys, and a number of full key pointers corresponding to the number offixed size partial keys, the method comprising: computing a search valuepartial key for the search value, wherein the search value partial keyis a fixed size and comprises a first byte offset and a subset of thesearch value; setting a current node to a root node and setting acurrent partial key to a first partial key of the root node, wherein thefirst partial key of the root node is the fixed size and comprises thefirst byte offset and a subset of a key referenced by a first full keypointer of the root node; and comparing the search value partial key tothe current partial key, wherein the comparing comprises; when thesearch value partial key is less than the current partial key, settingthe current node to a node referenced by a child group pointer of thecurrent node and setting the current partial key to a first partial keyof the identified node, wherein the first partial key of the identifiednode comprises a second byte offset and a subset of a key referenced bya first full key pointer of the identified node, and when the searchvalue partial key is greater than the current partial key, setting thecurrent partial key to a next partial key of the current node, whereinthe next partial key of the current node is the fixed size and comprisesa third byte offset and a subset of a key referenced by a next full keypointer of the current node, determining that the search value is notlocated in the database based on the comparing; and outputting anindication that the search value is not located in the database inresponse to the determining.